Performing a clear operation absent host intervention

ABSTRACT

Optimizations are provided for frame management operations, including a clear operation and/or a set storage key operation, requested by pageable guests. The operations are performed, absent host intervention, on frames not resident in host memory. The operations may be specified in an instruction issued by the pageable guests.

CROSS-REFERENCE TO RELATED APPLICATIONS

This application is a continuation of U.S. Ser. No. 14/576,729, entitled“EXECUTION OF A PERFORM FRAME MANAGEMENT FUNCTION INSTRUCTION,” filedDec. 19, 2014, which is a continuation of U.S. Ser. No. 14/136,086,entitled “EXECUTION OF A PERFORM FRAME MANAGEMENT FUNCTION INSTRUCTION,”filed Dec. 20, 2013 (U.S. Pat. No. 8,935,504, issued Jan. 13, 2015),which is a continuation of U.S. Ser. No. 13/941,887, entitled “EXECUTIONOF A PERFORM FRAME MANAGEMENT FUNCTION INSTRUCTION,” filed Jul. 15, 2013(U.S. Pat. No. 8,707,000, issued Apr. 22, 2014), which is a continuationof U.S. Ser. No. 13/554,056, entitled “EXECUTION OF A PERFORM FRAMEMANAGEMENT FUNCTION INSTRUCTION,” filed Jul. 20, 2012 (U.S. Pat. No.8,495,326, issued Jul. 23, 2013), which is a continuation of U.S. Ser.No. 13/292,160, entitled “CLEARING GUEST FRAMES ABSENT PAGING-IN TO HOSTMAIN STORAGE,” filed Nov. 9, 2011 (U.S. Pat. No. 8,239,649, issued Aug.7, 2012), which is a continuation of U.S. Ser. No. 12/036,725, entitled“OPTIMIZATIONS OF A PERFORM FRAME MANAGEMENT FUNCTION ISSUED BY PAGEABLEGUESTS,” filed Feb. 25, 2008 (U.S. Pat. No. 8,086,811, issued Dec. 27,2011), each of which is hereby incorporated herein by reference in itsentirety.

BACKGROUND

This invention relates, in general, to computing environments thatsupport pageable guests, and in particular, to facilitating processingof frame management operations within such environments.

Improving processing efficiency of computing environments continues tobe an important goal. One area in which improvements have been made, butfurther enhancements are needed, is in the area of supporting pageableguests.

An operating system frequently performs certain main storage (memory)management actions in support of its applications and to ensure securityand data integrity. These actions include clearing frames in real memoryand setting associated storage protection keys before the frames areassigned to an application or reassigned from one application toanother. When this operating system runs as a pageable guest, that is,in a virtual machine whose memory is paged by a hypervisor or host,these guest memory management operations may incur inefficiencies.

For example, if the host has paged out the contents of a guest framebeing reassigned, the clearing operation by the guest operating systemwould result in a host page fault, causing a context switch to the host.The host would then read the guest frame contents from host auxiliarystorage, and switch context back to the guest, which would then clearthe frame, immediately overlaying the contents which the host just readin.

Similarly, if the guest operating system clears a frame which anapplication has released and places it into an available pool to be usedlater, the resultant host page fault and handling will cause that guestframe to be made host resident, consuming host real memory at a timewhen no productive use will be made of that memory. The operation ofsetting the storage key on the guest frame commonly occurs at the sametime as clearing in these cases, and its handling entails manipulationof the same host interlock and data structures as the clearingoperation. Thus, treating frame management operations like clearing andkey setting the same as ordinary instructions leads to additionaloverhead for context switching between guest and host and for executingthe host page fault handler; wasted paging I/O bandwidth; less efficientuse of host memory; and repeated serialization and access to hosttranslation and control structures.

BRIEF SUMMARY

Thus, a need exists for a capability that reduces context switching andfacilitates processing in an environment that supports pageable guests.In one example, a need exists for a capability that providesoptimizations for frame management functions requested by pageableguests. For instance, a need exists for a capability that providesoptimizations for a memory clearing function (e.g., setting framecontents to zeros) issued by pageable guests. In a further example, aneed exists for a capability that provides optimizations for a set keyfunction issued by pageable guests. In one particular example, a needexists for a capability that provides optimizations when executing aPerform Frame Management Function instruction issued by a pageableguest, which could entail clearing guest frame contents and/or settingthe associated storage key or keys.

The shortcomings of the prior art are overcome and additional advantagesare provided through the provision of a computer system to facilitatemanagement of frames within a computing environment. The computer systemincludes a memory; and a processor in communication with the memory,wherein the computer system is configured to perform a method. Themethod includes based on initiating execution of an instruction issuedby a pageable guest executing within the computing environment andspecifying a clear operation, determining a usage intent characteristicfor a guest frame to be cleared, the guest frame not resident in hostmemory; and performing the clear operation absent host intervention, theperforming being based on the determined usage intent characteristic.

Methods and computer program products relating to one or more aspects ofthe present invention are also described and may be claimed herein.

Additional features and advantages are realized through the techniquesof the present invention. Other embodiments and aspects of the inventionare described in detail herein and are considered a part of the claimedinvention.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

One or more aspects of the present invention are particularly pointedout and distinctly claimed as examples in the claims at the conclusionof the specification. The foregoing and other objects, features, andadvantages of the invention are apparent from the following detaileddescription taken in conjunction with the accompanying drawings inwhich:

FIG. 1 depicts one embodiment of a computing environment to incorporateand use one or more aspects of the present invention;

FIG. 2 depicts one embodiment of an emulated computing environment toincorporate and use one or more aspects of the present invention;

FIG. 3 depicts one example of a frame descriptor used in accordance withan aspect of the present invention;

FIGS. 4A-4C depict examples of region table entries used in accordancewith an aspect of the present invention;

FIGS. 5A-5B depict examples of segment table entries used in accordancewith an aspect of the present invention;

FIG. 6 depicts one example of a page table entry used in accordance withan aspect of the present invention;

FIG. 7 depicts one example of a page status table entry used inaccordance with an aspect of the present invention;

FIG. 8A depicts one example of a format of a Perform Frame ManagementFunction (PFMF) instruction, in accordance with an aspect of the presentinvention;

FIG. 8B depicts one example of fields of registers specified in theinstruction of FIG. 8A, in accordance with an aspect of the presentinvention;

FIGS. 9A-9B depict one embodiment of the logic associated with thePerform Frame Management Function instruction, in accordance with anaspect of the present invention;

FIG. 10 depicts one embodiment of the logic to clear a frame for apageable guest, in accordance with an aspect of the present invention;

FIGS. 11A-11E depict one embodiment of the logic associated withperforming the Perform Frame Management Function instruction andoptimizations associated therewith, in accordance with an aspect of thepresent invention;

FIG. 12 depicts one example of various error processing associated withdynamic address translation, in accordance with an aspect of the presentinvention;

FIGS. 13A-13B depict one embodiment of the logic associated withperforming a resolve function, in accordance with an aspect of thepresent invention;

FIG. 14 depicts one embodiment of the logic associated with appending toa backing reclaim log, in accordance with an aspect of the presentinvention; and

FIG. 15 depicts one embodiment of a computer program productincorporating one or more aspects of the present invention.

DETAILED DESCRIPTION

In accordance with an aspect of the present invention, optimizations areprovided for frame management functions issued by pageable guests. Forexample, a capability is provided that enables a frame clearing functionin a pageable guest environment to be optimized. As another example, acapability is provided that enables a set key function issued bypageable guests to be optimized. As one particular example,optimizations are provided for a Perform Frame Management Function(PFMF) issued by pageable guests.

One embodiment of a computing environment to incorporate and use one ormore aspects of the present invention is described with reference toFIG. 1. Computing environment 100 is based, for instance, on thez/Architecture® offered by International Business Machines Corporation,Armonk, N.Y. The z/Architecture® is described in an IBM® publicationentitled, “z/Architecture Principles of Operation,” IBM® Publication No.SA22-7832-05, April, 2007, which is hereby incorporated herein byreference in its entirety. In one example, a computing environment basedon the z/Architecture® includes an eServer zSeries®, offered byInternational Business Machines Corporation, Armonk, N.Y. IBM®,z/Architecture® and zSeries® are registered trademarks of InternationalBusiness Machines Corporation, Armonk, N.Y., USA. Other names usedherein may be registered trademarks, trademarks, or product names ofInternational Business Machines Corporation or other companies.

As one example, computing environment 100 includes a central processorcomplex (CPC) 102 providing virtual machine support. CPC 102 includes,for instance, one or more virtual machines 104, one or more centralprocessors 106, at least one host 108 (e.g., a control program, such asa hypervisor), and an input/output subsystem 110, each of which isdescribed below. In this example, the virtual machines and host areincluded in memory.

The virtual machine support of the CPC provides the ability to operatelarge numbers of virtual machines, each capable of hosting a guestoperating system 112, such as Linux. Each virtual machine 104 is capableof functioning as a separate system. That is, each virtual machine canbe independently reset, host a guest operating system, and operate withdifferent programs. An operating system or application program runningin a virtual machine appears to have access to a full and completesystem, but in reality, only a portion of it is available.

In this particular example, the model of virtual machines is a V=Vmodel, in which the absolute or real memory of a virtual machine isbacked by host virtual memory, instead of real or absolute memory. Eachvirtual machine has a virtual linear memory space. The physicalresources are owned by host 108, and the shared physical resources aredispatched by the host to the guest operating systems, as needed, tomeet their processing demands. This V=V virtual machine (i.e., pageableguest) model assumes that the interactions between the guest operatingsystems and the physical shared machine resources are controlled by thehost, since the large number of guests typically precludes the host fromsimply partitioning and assigning the hardware resources to theconfigured guests. One or more aspects of a V=V model are furtherdescribed in an IBM® publication entitled “z/VM: Running Guest OperatingSystems,” IBM® Publication No. SC24-5997-02, October 2001, which ishereby incorporated herein by reference in its entirety.

Central processors 106 are physical processor resources that areassignable to a virtual machine. For instance, virtual machine 104includes one or more logical processors, each of which represents all ora share of a physical processor resource 106 that may be dynamicallyallocated to the virtual machine. Virtual machines 104 are managed byhost 108. As examples, the host may be implemented in microcode runningon processors 106 or be part of a host operating system executing on themachine. In one example, host 108 is a VM hypervisor, such as z/VM®,offered by International Business Machines Corporation, Armonk, N.Y. Oneembodiment of z/VM® is described in an IBM® publication entitled “z/VM:General Information Manual,” IBM Publication No. GC24-5991-04, October2001, which is hereby incorporated herein by reference in its entirety.

Input/output subsystem 110 directs the flow of information betweendevices and main storage. It is coupled to the central processingcomplex, in that it can be part of the central processing complex orseparate therefrom. The I/O subsystem relieves the central processors ofthe task of communicating directly with the I/O devices coupled to theCPC and permits data processing to proceed concurrently with I/Oprocessing.

In one embodiment, the host (e.g., z/VM®) and processor (e.g., System z)hardware/firmware interact with each other in a controlled cooperativemanner in order to process V=V guest operating system operations withoutrequiring transfer of control from/to the guest operating system and thehost. Guest operations can be executed directly without hostintervention via a facility that allows instructions to beinterpretively executed for a pageable storage mode guest. This facilityprovides an instruction, Start Interpretive Execution (SIE), which thehost can issue, designating a control block called a state descriptionwhich holds guest (virtual machine) state and controls. The instructionplaces the machine into an interpretive-execution mode in which guestinstructions and interruptions are processed directly, until a conditionrequiring host attention arises. When such a condition occurs,interpretive execution is ended, and either a host interruption ispresented, or the SIE instruction completes storing details of thecondition encountered; this latter action is called interception. Oneexample of interpretive execution is described in System/370 ExtendedArchitecture/Interpretive Execution, IBM Publication No. SA22-7095-01,September 1985, which is hereby incorporated herein by reference in itsentirety.

Another example of a computing environment to incorporate one or moreaspects of the present invention is depicted in FIG. 2. In this example,an emulated host computer system 200 is provided that emulates a hostcomputer 202 of a host architecture. In emulated host computer system200, a host processor (CPU) 204 is an emulated host processor (orvirtual host processor) and is realized through an emulation processor206 having a different native instruction set architecture than used bythe processors of host computer 202. Emulated host computer system 200has memory 208 accessible to emulation processor 206. In the exampleembodiment, memory 208 is partitioned into a host computer memory 210portion and an emulation routines 212 portion. Host computer memory 210is available to programs of emulated host computer 202 according to hostcomputer architecture, and may include both a host or hypervisor 214 andone or more virtual machines 216 running guest operating systems 218,analogous to the like-named elements in FIG. 1.

Emulation processor 206 executes native instructions of an architectedinstruction set of an architecture other than that of the emulatedprocessor 204. The native instructions are obtained, for example, fromemulation routines memory 212. Emulation processor 206 may access a hostinstruction for execution from a program in host computer memory 210 byemploying one or more instruction(s) obtained in a sequence &access/decode routine which may decode the host instruction(s) accessedto determine a native instruction execution routine for emulating thefunction of the host instruction accessed. One such host instruction maybe, for example, a Start Interpretive Execution (SIE) instruction, bywhich the host seeks to execute a guest program in a virtual machine.The emulation routines 212 may include support for this instruction, andfor executing a sequence of guest instructions in a virtual machine 216in accordance with the definition of this SIE instruction.

Other facilities that are defined for host computer system 202architecture may be emulated by architected facilities routines,including such facilities as general purpose registers, controlregisters, dynamic address translation, and I/O subsystem support andprocessor cache for example. The emulation routines may also takeadvantage of functions available in emulation processor 206 (such asgeneral registers and dynamic translation of virtual addresses) toimprove performance of the emulation routines. Special hardware andoffload engines may also be provided to assist processor 206 inemulating the function of host computer 202.

In providing the optimizations described herein, reference is made tovarious facilities and data structures (e.g., tables, lists). Examplesof these facilities and data structures are described below tofacilitate an understanding of one or more aspects of the presentinvention.

One facility that is referenced is dynamic address translation (DAT) andenhanced dynamic address translation (EDat). Dynamic address translationis the process of translating a virtual address during a storagereference into the corresponding real address or absolute address. Thevirtual address may be a primary virtual address, a secondary virtualaddress, an Access Register specified virtual address, or a home virtualaddress. These addresses are translated by means of a primary, asecondary, an AR specified, or a home address space control element,respectively. After selection of the appropriate address space controlelement, the translation process is the same for all of the four typesof virtual address. DAT may use from five to two levels of tables(region first table, region second table, region third table, segmenttable, and page table) as transformation parameters. Enhanced DAT mayuse from five to one levels of table, by omitting the page table forsome or all translations. The designation (origin and length) of thehighest level table for a specific address space is called an addressspace control element, and it is found for use by DAT in a controlregister or as specified by an access register. Alternatively, theaddress space control element for an address space may be a real spacedesignation, which indicates that DAT is to translate the virtualaddress simply by treating it as a real address and without using anytables.

DAT uses, at different times, the address space control elements indifferent control registers or specified by the access registers. Thechoice is determined by the program-specified translation mode specifiedin the current PSW (Program Status Word). Four translation modes areavailable: primary space mode, secondary space mode, access registermode, and home space mode. Different address spaces are addressabledepending on the translation mode.

The result of enhanced DAT upon a virtual address may be either a realor an absolute address. If it is a real address, a prefixing operationis then performed to obtain the corresponding absolute address, whichcan be used to reference memory. Prefixing provides the ability toassign the range of real addresses 0-8191 (as an example) to a differentarea in absolute storage for each CPU, thus permitting more than one CPUsharing main storage to operate concurrently with a minimum ofinterference, especially in the processing of interruptions. Prefixingcauses real addresses in the range 0-8191 to correspond one-for-one tothe area of 8K byte absolute addresses (the prefix area) identified bythe value in bit positions 0-50 of the prefix register for the CPU, andthe area of real addresses identified by that value in the prefixregister to correspond one-for-one to absolute addresses 0-8191. Theremaining real addresses are the same as the corresponding absoluteaddresses. This transformation allows each CPU to access all of mainstorage, including the first 8K bytes and the locations designated bythe prefix registers of other CPUs.

Dynamic address translation, prefixing, and enhanced DAT are describedin more detail in U.S. Publication No. 2009/0187724A1, entitled,“Dynamic Address Translation with Frame Management,” Greiner et al.,(IBM Docket No.: POU920070313US1), published Jul. 23, 2009, which ishereby incorporated herein by reference in its entirety.

Another facility referenced is a Host Page Management Assist (HPMA)facility that includes various functions that can be invoked during theinterpretation of a pageable storage mode guest. One function that isinvoked is a resolve host page function used to dynamically resolve ahost page invalid condition. One example of HPMA is described in U.S.Publication No. 2005/0268071A1, entitled “Facilitating Management ofStorage of a Pageable Mode Virtual Environment Absent Intervention of aHost of the Environment,” Blandy et al., published Dec. 1, 2005, whichis hereby incorporated herein by reference in its entirety.

A yet further facility referenced is a Collaborative Memory Management(CMM) facility that provides a vehicle for communicating granular pagestate information between a pageable guest and its host. It includes abacking reclaim log (CBRL) used to hold a list of frames backing unusedpages. One example of CMM is described in U.S. Patent ApplicationPublication No. US 2007/0016904 A1, entitled, “Facilitating ProcessingWithin Computing Environment Supporting Pageable Guests,” Adlung et al.,published Jan. 18, 2007, which is hereby incorporated herein byreference in its entirety.

Reference is also made to various structures, which are described below.

Frame Descriptor

A frame descriptor describes a host page frame; that is, an area of realmemory (frame) capable of holding a portion of virtual memory (page). Itis allocated, deallocated, and initialized by the host and may beupdated by Host Page Management Assist functions (as described, forinstance in U.S. Publication No. 2005/0268071A1, entitled “FacilitatingManagement of Storage of a Pageable Mode Virtual Environment AbsentIntervention of a Host of the Environment,” Blandy et al., publishedDec. 1, 2005, which is hereby incorporated herein by reference in itsentirety).

In one example, a frame descriptor 300 (FIG. 3) is, for instance, a32-byte block residing in host home space virtual storage on a 32 byteboundary, and includes the following fields, as examples:

-   -   (a) Next Frame Descriptor Address 302: In one example, the        contents of this field, with five zeros appended on the right,        specify the host home space virtual address of the next frame        descriptor on the list. A value of zero indicates that the frame        descriptor is the last on the list.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.    -   (b) Page Frame Real Address or PTE Copy 304: When the frame        descriptor is in the available frame descriptor list (AFDL), the        contents of this field, with twelve zeros appended on the right,        specify the host real address of the first byte (byte 0) of a        host frame that is available for allocation to provide host        storage.    -   When the frame descriptor is in a processed frame descriptor        list (PFDL), this field includes a copy of the page table entry        (PTE) designated by the page table entry address field, as it        appeared before the host page was resolved.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.    -   (c) Page Table Entry Address 306: When the frame descriptor is        on the processed frame descriptor list, the contents of this        field, with three zeros appended on the right, specify the host        real or host absolute address of the page table entry for the        host virtual page.    -   This field is initialized by the host and may be changed by the        host or by Host Page Management Assist functions.

Multiple frame descriptors may be linked to one another to form a list,such as an available frame descriptor list (AFDL) or a processed framedescriptor list (PFDL). A frame descriptor exists in one of the twolists. A separate pair of these lists is provided for each CPU. Theorigins of the AFDL and PFDL for a CPU are designated by means of fieldsin the prefix area of the CPU.

The available frame descriptor list (AFDL) is a list of framedescriptors that describes host frames the host has cleared and has madeavailable for allocation to host pages. The AFDL is designated by anAFDL origin (AFDLO) at a specified host real address.

The contents of the AFDLO, with five zeros appended on the right,specify the host home space virtual address of the first framedescriptor on the AFDL. A value of zero indicates that the list isempty.

The AFDLO is initialized by the host and may be changed by the host orHost Page Management Assist functions. The AFDLO is changed, in oneembodiment, by means of a non-interlocked update operation.

The processed frame descriptor list (PFDL) is a list of framedescriptors that describes host frames that have been used to resolvehost page invalid conditions during guest interpretation. The hostframes that are described by the PFDL have been assigned to host pagesthat provide storage for a guest. The PFDL is designated by a PFDLorigin (PFDLO) at a specified host real address. The contents of thePFDLO, with five zeros appended on the right, specify the host homespace virtual address of the first frame descriptor on the PFDL. A valueof zero indicates that the list is empty.

The PFDLO is initialized by the host and may be changed by the host or aHost Page Management Assist function. The PFDLO is changed, in oneembodiment, by means of a doubleword concurrent interlocked updateoperation that maintains the integrity of the list.

Region Table Entries

The term “region table entry” indicates a region first table entry, aregion second table entry, or a region third table entry. The level(first, second, or third) of the table containing an entry is identifiedby the table type (TT) bits in the entry. Examples of the formats ofentries fetched from the region first table, region second table, andregion third table are depicted in FIGS. 4A-4C. In particular, FIG. 4Adepicts one embodiment of the format of a Region First Table entry 400;FIG. 4B depicts one embodiment of the format of a Region Second Tableentry 430; and FIG. 4C depicts one embodiment of the format of a RegionThird Table entry 460.

As examples, the fields in the three levels of region table entries areallocated as follows:

Region Second Table Origin 402, Region Third Table Origin 432, andSegment Table Origin 462: A region first table entry includes a regionsecond table origin. A region second table entry includes a region thirdtable origin. A region third table entry includes a segment tableorigin. The following description applies to each of the three origins.In one example, bits 0-51 of the entry, with 12 zeros appended on theright, form a 64-bit address that designates the beginning of the nextlower level table.

DAT Protection Bit (P) 406, 436, 466: When enhanced DAT applies, bit 54is treated as being OR'ed with the DAT protection bit in each subsequentregion table entry, segment table entry, and, when applicable, pagetable entry used in the translation. Thus, when the bit is one, DATprotection applies to the entire region or regions specified by theregion table entry. When the enhanced DAT facility is not installed, orwhen the facility is installed but the enhanced DAT enablement controlis zero, bit 54 of the region table entry is ignored.

Region Second Table Offset 408, Region Third Table Offset 438, andSegment Table Offset (TF) 468: A region first table entry includes aregion second table offset. A region second table entry includes aregion third table offset. A region third table entry includes a segmenttable offset. The following description applies to each of the threeoffsets. Bits 56 and 57 of the entry specify the length of a portion ofthe next lower level table that is missing at the beginning of thetable; that is, the bits specify the location of the first entryactually existing in the next lower level table. The bits specify thelength of the missing portion in units of 4,096 bytes, thus making thelength of the missing portion variable in multiples of 512 entries. Thelength of the missing portion, in units of 4,096 bytes, is equal to theTF value. The contents of the offset field, in conjunction with thelength field, bits 62 and 63, are used to establish whether the portionof the virtual address (RSX, RTX, or SX) to be translated by means ofthe next lower level table designates an entry that actually exists inthe table.

Region Invalid Bit (I) 410, 440, 470: A region is a contiguous range of,for example, 2 gigabytes of virtual addresses. Bit 58 in a region firsttable entry or region second table entry controls whether the set ofregions associated with the entry is available. Bit 58 in a region thirdtable entry controls whether the single region associated with the entryis available. When bit 58 is zero, address translation proceeds by usingthe region table entry. When the bit is one, the entry cannot be usedfor translation.

Table Type Bits (TT) 412, 442, 472: Bits 60 and 61 of the region firsttable entry, region second table entry, and region third table entryidentify the level of the table containing the entry, as follows: Bits60 and 61 identify the correct table level, considering the type oftable designation that is the address space control element being usedin the translation and the number of table levels that have so far beenused; otherwise, a translation specification exception is recognized. Asan example, the following table shows the table type bits:

Table Type bits for region table Entries Bits 60 and 61 Region-TableLevel 11 First 10 Second 01 Third

Region Second Table Length 414, Region Third Table Length 444, andSegment Table Length 474 (TL): A region first table entry includes aregion second table length. A region second table entry includes aregion third table length. A region third table entry includes a segmenttable length. The following description applies to each of the threelengths. Bits 62 and 63 of the entry specify the length of the nextlower level table in units of 4,096 bytes, thus making the length of thetable variable in multiples of 512 entries. The length of the next lowerlevel table, in units of 4,096 bytes, is one more than the TL value. Thecontents of the length field, in conjunction with the offset field, bits56 and 57, are used to establish whether the portion of the virtualaddress (RSX, RTX, or SX) to be translated by means of the next lowerlevel table designates an entry that actually exists in the table.

All other bit positions of the region table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the region table entry shouldcontain zeros even if the table entry is invalid.

Segment Table Entries

When enhanced DAT does not apply, or when enhanced DAT applies and theSTE format control, bit 53 of the segment table entry is zero, the entryfetched from the segment table has the format (e.g., Format 0) asdepicted in FIG. 5A. When enhanced DAT applies and the STE formatcontrol is one, the entry fetched from the segment table has, forexample, the format (e.g., Format 1) as depicted in FIG. 5B.

As one example, a Format 0 segment table entry 500 (FIG. 5A) includesthe following fields:

Page Table Origin 502: When enhanced DAT does not apply, or whenenhanced DAT applies but the STE format control, bit 53 of the segmenttable entry, is zero, bits 0-52, with 11 zeros appended on the right,form a 64-bit address that designates the beginning of a page table. Itis unpredictable whether the address is real or absolute.

STE Format Control (FC) 506: When enhanced DAT applies, bit 53 is theformat control for the segment table entry, as follows:

-   -   When the FC bit is zero, bits 0-52 of the entry form the page        table origin, and bit 55 is reserved.    -   When the FC bit is one, bits 0-43 of the entry form the segment        frame absolute address, bit 47 is the ACCF validity control,        bits 48-51 are the access control bits, bit 52 is the fetch        protection bit, and bit 55 is the change recording override.        When enhanced DAT does not apply, bit 53 is ignored.

DAT Protection Bit (P) 508: Bit 54, when one, indicates that DATprotection applies to the entire segment.

-   -   When enhanced DAT does not apply, bit 54 is treated as being        OR'ed with the DAT protection bit in the page table entry used        in the translation.    -   When enhanced DAT applies, the DAT protection bit in any and all        region table entries used in the translation are treated as        being OR'ed with the DAT protection bit in the segment table        entry; when the STE format control is zero, the DAT protection        bit in the STE is further treated as being OR'ed with the DAT        protection bit in the page table entry.

Segment Invalid Bit (I) 510: Bit 58 controls whether the segmentassociated with the segment table entry is available.

-   -   When the bit is zero, address translation proceeds by using the        segment table entry.    -   When the bit is one, the segment table entry cannot be used for        translation.

Common Segment Bit (C) 512: Bit 59 controls the use of the translationlookaside buffer (TLB) copies of the segment table entry. When enhancedDAT does not apply or when enhanced DAT applies but the format controlis zero, bit 59 also controls the use of the TLB copies of the pagetable designated by the segment table entry.

-   -   A zero identifies a private segment; in this case, the segment        table entry and any page table it designates may be used only in        association with the segment table origin that designates the        segment table in which the segment table entry resides.    -   A one identifies a common segment; in this case, the segment        table entry and any page table it designates may continue to be        used for translating addresses corresponding to the segment        index, even though a different segment table is specified.

However, TLB copies of the segment table entry and any page table for acommon segment are not usable if the private space control, bit 55, isone in the address space control element used in the translation or ifthat address space control element is a real space designation. Thecommon segment bit is to be zero if the segment table entry is fetchedfrom storage during a translation when the private space control is onein the address space control element being used; otherwise, atranslation specification exception is recognized.

Table Type Bits (TT) 514: Bits 60 and 61 of the segment table entry are00 binary to identify the level of the table containing the entry. Themeanings of possible values of bits 60 and 61 in a region table entry orsegment table entry are as follows:

Table Type Bits 60, 61 Bits 60 and 61 Table Level 11 Region-first 10Region-second 01 Region-third 00 Segment

Bits 60 and 61 are to identify the correct table level, considering thetype of table designation that is the address space control elementbeing used in the translation and the number of table levels that haveso far been used; otherwise, a translation specification exception isrecognized.

All other bit positions of the segment table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the segment table entry shouldcontain zeros even if the table entry is invalid.

As one example, a Format 1 segment table entry 550 (FIG. 5B) includesthe following fields:

Segment Frame Absolute Address (SFAA) 552: When enhanced DAT applies andthe STE format control is one, bits 0-43 of the entry, with 20 zerosappended on the right, form the 64-bit absolute address of the segment.

ACCF Validity Control (AV) 556: When enhanced DAT applies and the STEformat control is one, bit 47 is the access control bits and fetchprotection bit (ACCF) validity control. When the AV control is zero,bits 48-52 of the segment table entry are ignored. When the AV controlis one, bits 48-52 are used as described below.

Access Control Bits (ACC) 558: When enhanced DAT applies, the STE formatcontrol is one, and the AV control is one, bits 48-51 of the segmenttable entry include the access control bits that may be used for any keycontrolled access checking that applies to the address.

Fetch Protection Bit (F) 560: When enhanced DAT applies, the STE formatcontrol is one, and the AV control is one, bit 52 of the segment tableentry includes the fetch protection bit that may be used for any keycontrolled access checking that applies to the address.

STE Format Control (FC) 562: When enhanced DAT applies, bit 53 is theformat control for the segment table entry, as follows:

-   -   When the FC bit is zero, bits 0-52 of the entry form the page        table origin, and bit 55 is reserved.    -   When the FC bit is one, bits 0-43 of the entry form the segment        frame absolute address, bit 47 is the ACCF validity control,        bits 48-51 are the access control bits, bit 52 is the fetch        protection bit, and bit 55 is the change recording override.        When enhanced DAT does not apply, bit 53 is ignored.

DAT Protection Bit (P) 564: Bit 54, when one, indicates that DATprotection applies to the entire segment.

-   -   When enhanced DAT does not apply, bit 54 is treated as being        OR'ed with the DAT protection bit in the page table entry used        in the translation.    -   When enhanced DAT applies, the DAT protection bit in any and all        region table entries used in the translation are treated as        being OR'ed with the DAT protection bit in the segment table        entry; when the STE format control is zero, the DAT protection        bit in the STE is further treated as being OR'ed with the DAT        protection bit in the page table entry.

Change Recording Override (CO) 566: When enhanced DAT applies, and theSTE format control is one, bit 55 of the segment table entry is thechange recording override for the segment. When enhanced DAT does notapply, or when enhanced DAT applies but the STE format control is zero,bit 55 of the segment table entry is ignored.

Segment Invalid Bit (I) 568: Bit 58 controls whether the segmentassociated with the segment table entry is available.

-   -   When the bit is zero, address translation proceeds by using the        segment table entry.    -   When the bit is one, the segment table entry cannot be used for        translation.

Common Segment Bit (C) 570: Bit 59 controls the use of the translationlookaside buffer (TLB) copies of the segment table entry. When enhancedDAT does not apply or when enhanced DAT applies but the format controlis zero, bit 59 also controls the use of the TLB copies of the pagetable designated by the segment table entry.

-   -   A zero identifies a private segment; in this case, the segment        table entry and any page table it designates may be used only in        association with the segment table origin that designates the        segment table in which the segment table entry resides.    -   A one identifies a common segment; in this case, the segment        table entry and any page table it designates may continue to be        used for translating addresses corresponding to the segment        index, even though a different segment table is specified.

However, TLB copies of the segment table entry and any page table for acommon segment are not usable if the private space control, bit 55, isone in the address space control element used in the translation or ifthat address space control element is a real space designation. Thecommon segment bit is to be zero if the segment table entry is fetchedfrom storage during a translation when the private space control is onein the address space control element being used; otherwise, atranslation specification exception is recognized.

Table Type Bits (TT) 572: Bits 60 and 61 of the segment table entry are00 binary to identify the level of the table containing the entry. Themeanings of possible values of bits 60 and 61 in a region table entry orsegment table entry are as follows:

Table Type Bits 60, 61 Bits 60 and 61 Table Level 11 Region-first 10Region-second 01 Region-third 00 Segment

Bits 60 and 61 are to identify the correct table level, considering thetype of table designation that is the address space control elementbeing used in the translation and the number of table levels that haveso far been used; otherwise, a translation specification exception isrecognized.

All other bit positions of the segment table entry are reserved forpossible future extensions and should contain zeros; otherwise, theprogram may not operate compatibly in the future. When enhanced DATapplies, the reserved bit positions of the segment table entry shouldcontain zeros even if the table entry is invalid.

Page Table Entries

The state information for guest blocks (e.g., an area (e.g., 4 K-Bytes)in absolute memory that has associated therewith a single storage keyand CMM state) is maintained, for instance, in host page tables (PTs)and page status tables (PGSTs) that describe a guest's memory. Thesetables include, for instance, one or more page table entries (PTEs) andone or more page status table entries (PGSTEs), respectively, which aredescribed in further detail below.

One example of a page table entry 600 is described with reference toFIG. 6. In one embodiment, the fields in the page table entry areallocated as follows:

Page Frame Real Address (PFRA) 602: Bits 0-51 provide the leftmost bitsof a real (in this case host real) storage address. When these bits areconcatenated with the 12-bit byte index field of the virtual address onthe right, a 64-bit real address is obtained.

Page Invalid Bit (I) 604: Bit 53 controls whether the page associatedwith the page table entry is available. When the bit is zero, addresstranslation proceeds by using the page table entry. Further, with regardto collaborative memory management (CMM) between host and guest, thehost state is r (resident; i.e., the guest block is present in a hostframe). When the bit is one, the page table entry is not used fortranslation, and the CMM host state is p (preserved; i.e., the guestblock is not present in a host frame, but has been preserved by the hostin some auxiliary storage) or z (logically zero; i.e., the guest blockis not present in a host frame and the contents of the guest block areknown to be zeros), as determined by PGSTE.Z.

DAT Protection Bit (P) 606: Bit 54 controls whether store accesses canbe made in the page. This protection mechanism is in addition to the keycontrolled protection and low address protection mechanisms. The bit hasno effect on fetch accesses. If the bit is zero, stores are permitted tothe page, subject to the following additional constraints:

-   -   The DAT protection bit being zero in the segment table entry        used in the translation.    -   When enhanced DAT applies, the DAT protection bit being zero in        all region table entries used in the translation.

If the DAT protection bit is one, stores are disallowed. When no higherpriority exception conditions exist, an attempt to store when the DATprotection bit is one causes a protection exception to be recognized.The DAT protection bit in the segment table entry is treated as beingOR'ed with bit 54 when determining whether DAT protection applies to thepage. When enhanced DAT applies, the DAT protection bit in any regiontable entries used in translation are also treated as being OR'ed withbit 54 when determining whether DAT protection applies.

Other protection mechanisms, such as key-controlled protection,low-address protection, and access-list-controlled protection, may applyindependently of DAT protection and may also prohibit accesses.

Change Recording Override (CO) 608: When enhanced DAT does not apply,bit 55 of the page table entry is to contain zero; otherwise, atranslation specification exception is recognized as part of theexecution of an instruction using that entry for address translation.When enhanced DAT applies and the STE format control is zero, bit 55 ofthe page table entry is the change recording override for the page.

In addition to the above, in one example, bit position 52 of the entryis to contain zero; otherwise, a translation specification exception isrecognized as part of the execution of an instruction using that entryfor address translation. Bit positions 56-63 are not assigned and areignored.

One example of a page status table entry is described with reference toFIG. 7. A page status table entry 700 includes, for instance, thefollowing:

-   -   (a) Acc 702: Access control key;    -   (b) FP 704: Fetch protection indicator;    -   (c) Page Control Interlock (PCL) 706: This is the interlock        control for serializing updates to a page table entry (PTE) and        corresponding PGSTE, except for the PGSTE status area and PGSTE        bits that are marked as reserved.    -   (d) HR 708: Host reference backup indicator;    -   (e) HC 710: Host change backup indicator;    -   (f) GR 712: Guest reference backup indicator;    -   (g) GC 714: Guest change backup indicator;    -   (h) Status 716: Intended for host program use.    -   (i) Page Content Logically Zero Indicator (Z) 718: This bit is        meaningful when the corresponding PTE page invalid (PTE.I) bit        is one.        -   When Z is one, the content of the page that is described by            this PGSTE and corresponding PTE is considered to be zero.            Any prior content of the page does not have to be preserved            and may be replaced by a page of zeros.        -   When Z is zero, the content of the page described by the            PGSTE and corresponding PTE is not considered to be zero.            The content of the page is preserved by the host.        -   When the Z bit is one and the corresponding PTE.I bit is            one, the CMM host state is z (logically zero). This means            that the page content may be replaced by the host or by a            function of the Host Page Management Assist facility.        -   When the Z bit is one, the corresponding PTE.I bit is one,            and the page content is replaced, the page should be            replaced by associating it with a frame that has been set to            zeros.        -   When the Z bit is zero and the PTE invalid bit is one, the            CMM host state is p (preserved).    -   (j) Page Class (PC) 720: When zero, the page described by the        PGSTE and corresponding PTE is a class 0 page and the delta        pinned page count array (DPPCA) for class 0 pages is used for        counting pinning and unpinning operations for the page. When        one, the page described by the PGSTE and corresponding PTE is a        class 1 page and the DPPCA for class 1 pages is used for        counting pinning and unpinning operations for the page.    -   (k) Pin Count Overflow (PCO) 722: When one, the pin count field        is in an overflow state. In this case, the total pin count is        kept by the host in another data structure not accessed by the        machine. When zero, the pin count field is not in an overflow        state.    -   (l) Frame Descriptor On Processed Frame Descriptor List (FPL)        724: When one, a frame descriptor for the page described by the        PGSTE and corresponding PTE is in a processed frame descriptor        list. The frame descriptor identifies the host frame that was        used by a HPMA resolve host page function for the page.    -   (m) Page Content Replacement Requested (PCR) 726: When one, page        content replacement was requested when the HPMA resolve host        page function was invoked for the page represented by the PGSTE        and corresponding PTE.    -   (n) Usage State (US) 728: For collaborative memory management        between host and guest, this field indicates whether the guest        state is S (stable; i.e., the contents of a stable block remain        equal to what was set by the guest); U (unused; i.e., the        contents of an unused block are not meaningful to the guest); V        (volatile; i.e., the contents of a volatile block are meaningful        to the guest, but the host may at any time discard the contents        of the block and reclaim the backing frame); or P (potentially        volatile; i.e., the contents of a potentially volatile block are        meaningful to the guest, but based upon guest change history,        the host either may discard or should preserve the contents of        the block).    -   (o) Status 730: Intended for host program use.    -   (p) Pin Count 732: An unsigned binary integer count used to        indicate whether the content of the host virtual page        represented by the PGSTE and corresponding PTE is pinned in the        real host frame specified by the page frame real address field        of the PTE. When the value of this field is greater than zero or        the page count overflow (PCO) bit is one, the corresponding page        is considered to be pinned. When the value of this field is zero        and the PCO bit is zero, the corresponding page is not        considered to be pinned.        -   At the time a page is pinned by either the host or the CPU,            this field should be incremented by 1. At the time a page is            unpinned by either the host or the CPU, this field should be            decremented by 1.        -   When the value of the pin count field is greater than zero            or the PCO bit is one, the corresponding PTE.I (page            invalid) bit is to be zero. Otherwise, unpredictable results            may occur.        -   While a page is pinned, the host program should not change            the contents of the PTE page frame real address (PFRA)            field, the setting of the PTE page invalid (I) bit, or the            setting of the page protection (P) bit in the PTE or segment            table entry (STE). Otherwise unpredictable results may            occur.

Further details regarding page table entries and page tables, as well assegment table entries, are provided in an IBM® publication entitled,“z/Architecture Principles of Operation,” IBM® Publication No.SA22-7832-05, April 2007, which is hereby incorporated herein byreference in its entirety. Moreover, further details regarding the PGSTEare described in U.S. Publication No. 2005/0268071A1, entitled“Facilitating Management of Storage of a Pageable Mode VirtualEnvironment Absent Intervention of a Host of the Environment,” Blandy etal., published Dec. 1, 2005; and in U.S. Patent Application PublicationNo. US 2007/0016904 A1, entitled, “Facilitating Processing WithinComputing Environment Supporting Pageable Guests,” Adlung et al.,published Jan. 18, 2007, each of which is hereby incorporated herein byreference in its entirety.

In one embodiment, there is one page status table per page table, thepage status table is the same size as the page table, a page statustable entry is the same size as a page table entry, and the page statustable is located at a fixed displacement (in host real memory) from thepage table. Thus, there is a one-to-one correspondence between each pagetable entry and page status table entry. Given the host's virtualaddress of a page, both the machine and the host can easily locate thepage status table entry that corresponds to a page table entry for aguest block.

Perform Frame Management Function (PFMF)

In accordance with an aspect of the present invention, optimizations areprovided for one or more aspects of a Perform Frame Management Function.Thus, prior to describing the optimizations, one example of the PFMFinstruction and processing associated therewith are described. This isto facilitate an understanding of one or more aspects of the presentinvention.

As one example, a format of a Perform Frame Management Function isdescribed with reference to FIG. 8A. As depicted, a PFMF instruction 800includes an opcode 802 identifying the PFMF instruction, a firstregister field 804 including a first operand, and a second registerfield 806 indicating a second operand address.

Subject to the controls in the first operand register, a framemanagement function is performed for the storage frame designated by thesecond operand address.

One example of the contents of general register R1 804 (R1 designates ageneral register) are described with reference to FIG. 8B and include,for instance:

Frame Management Function Indications 810: Bit positions 44-47 ofgeneral register R1 include the frame management function indications(FMFI), as follows:

-   -   Set Key Control (SK) 812: Bit 46 controls whether the storage        key for each 4K-byte block in the frame is set from bits 56-62        of general register R1. When the SK control is zero, the keys        are not set; when the SK control is one, the keys are set.    -   Clear Frame Control (CF) 814: Bit 47 controls whether the        contents of the frame are set to zeros. When the CF control is        zero, no clearing operation is performed. When the CF control is        one, the frame is cleared to zeros.

Usage Indication (UI) 816: Bit position 48 of general register R1includes the usage indication (UI). When bit 48 is zero, it indicatesthat the program does not anticipate immediate usage of the frame. Whenbit 48 is one, it indicates that the program anticipates usage of theframe in the near future. The definition of immediate or near future isa program dependent indication of likelihood of the program using theframe in the immediate or near future. That is, it is a program definedtime. For instance, if the guest operating system is using PFMF toprovide memory requested by an application, for instance through aGETMAIN or malloc( )request, then immediate usage is likely, so UI wouldbe set to one. If the application has returned the memory which theframe backs, for instance through a FREEMAN or free( )request, and theoperating system is using PFMF to clear the frame prior to placing itinto an available pool, imminent usage is not likely, so UI would be setto zero. Other examples are also possible.

Frame Size Code (FSC) 818: Bits 49-51 of general register R1 include theframe size code (FSC), as follows:

Meaning of frame size codes FSC Meaning 0 4K-byte frame 1 1M-byte frame2-7 Reserved

Reference Bit Update Mask (MR) 820: When the set key control, bit 46 ofgeneral register R1, is one, bit 53 of general register R1 controlswhether updates to the reference bit in the storage key may be bypassed.

Change Bit Update Mask (MC) 822: When the set key control, bit 46 ofgeneral register R1, is one, bit 54 of general register R1 controlswhether updates to the change bit in the storage key may be bypassed.

In one embodiment, the handling of the MR and MC bits is the same as thehandling of the corresponding bits of the M3 field of a Set Storage KeyExtended instruction, except that general register R1 is not updatedwith the contents of the previous key, and the condition code is notchanged.

Key 824: When the set key control, bit 46 of general register R1, isone, bits 56-62 of the register include the storage key to be set foreach 4K-byte block in the frame, with the access protection bits 826,fetch protection bit 828, reference bit 830, and change bit 832 in bitpositions 56-59, 60, 61, and 62, respectively.

General register R2 806, examples of which are also depicted in FIG. 8B,includes, for instance, the real or absolute address 850 of the storageframe upon which the frame management function is to be performed. Whenthe frame size code designates a 4K-byte block, the second operandaddress is real; when the frame size code designates a 1M-byte block thesecond operand address is absolute. The handling of the address ingeneral register R2 depends on the addressing mode. In the 24-bitaddressing mode, the contents of bit positions 40-51 of the register,with 12 rightmost zeros appended, are the address, and bits 0-39 and52-63 in the register are ignored. In the 31-bit addressing mode, thecontents of bit positions 33-51 of the register, with 12 rightmost zerosappended, are the address, and bits 0-32 and 52-63 in the register areignored. In the 64-bit addressing mode, the contents of bit positions0-51 of the register, with 12 rightmost zeros appended, are the address,and bits 52-63 in the register are ignored.

In processing a PFMF instruction, the following occurs, in one example:

When the frame size code is 0, the specified frame management functionsare performed for the 4K-byte frame specified by the second operand.General register R2 is unmodified in this case.

When the frame size code is 1, the specified frame management functionsare performed for one or more 4K-byte blocks within the 1M-byte frame,beginning with the block specified by the second operand address, andcontinuing to the right with each successive block up to the next1M-byte boundary. In this case, the Perform Frame Management Function isinterruptible, and processing is as follows:

-   -   When an interruption occurs (other than one that follows        termination), the second operand address in general register R2        is updated by the number of 4K-byte blocks processed, so the        instruction, when re-executed, resumes at the point of        interruption.    -   When the instruction completes without interruption, the second        operand address in general register R2 is updated to the next        1M-byte boundary.

When the frame size code is 1 in the 24-bit or 31-bit addressing modes,the leftmost bits which are not part of the address in bit positions32-63 of general register R2 are set to zeros; bits 0-31 of the registerare unchanged.

When the clear frame control is one, references to main storage withinthe second operand are not necessarily single access references and arenot necessarily performed in a left-to-right direction as observed byother CPUs and by channel programs. The clear operation is not subjectto key controlled protection.

When the storage key control is one, the operation for each 4K-byteblock is similar to that described in relation to the SSKE except thatwhen the keys for multiple blocks are set, the condition code and thecontents of general register R1 are unchanged.

A serialization and checkpoint synchronization function is performedbefore the operation begins and again after the operation is completed,except that when the seven bits of the storage keys to be set are thesame as bits 56-62 of general register R1, or when the MR and MC bitsallow the storage keys to remain unchanged, it is unpredictable whetherthe serialization and checkpoint synchronization operations areperformed after the operation completes. It is unpredictable whether theclear frame or the set key operation is performed first when both of therespective controls are one. Provided that there is no other access tothe storage by other CPUs or a channel subsystem, the final results ofthe instruction reflect the specified key value including the specifiedR and C values when MR and MC are zero.

Special Conditions

A specification exception is recognized and the operation is suppressedfor any of the following conditions:

-   -   Bits 32-45, 52, 55, or 63 of general register R1 are not zero.    -   The frame size code specifies a reserved value.

Condition Code: The code remains unchanged.

Perform Frame Management Function (PFMF)—Clear Frame

One embodiment of the logic of the Perform Frame Management Function inwhich the indicated frames are cleared is described with reference toFIG. 9A.

Generally, this embodiment involves, at 900, obtaining a machineinstruction defined for the machine architecture. The instructionincludes an opcode for a frame management instruction. The machineinstruction has a first field identifying a first general register and asecond field identifying a second general register. At 902, clear frameinformation having a frame size field is obtained from the first generalregister. At 904, a determination is made whether the frame size fieldindicates that a storage frame is one of a small block of data in memoryor a large block of data in memory. Then, at 906, a second operandaddress of the storage frame upon which the machine instruction is to beperformed is obtained from the second general register. The secondoperand address being either a real address of the small block of datain memory or an absolute address of the large block of data in memory.At 908, if the indicated storage frame is a small block of data, thenthe small block of data is cleared by setting the bytes of the smallblock of data to zero. At 910, if the indicated storage frame is a largeblock of data, an operand address of an initial first block of data ofthe large block of data is obtained from the second general register.The large block of data having a first plurality of first blocks ofdata. A second plurality of the first plurality of first blocks of datais cleared by setting the bytes of the second plurality of the firstplurality of first blocks of data to zero.

The embodiment further involving obtaining a block size indicator from afield of the machine instruction. Based on the block size indicator, adetermination is made whether the addressed operand is either the largeblock of data or the small block of data. The small block of data beinga same size as the first block of data. For the large block of data, anaddress of a next block of data is saved in the second general register.The next block of data being a block of data following the firstplurality of first blocks of data. For the large block of data, theaddress of the next block of data being determined by any one of a)encountering a boundary of the large block of data or b) responding to aprogram interruption event.

In another embodiment, clear frame information is obtained from thefirst general register. The clear frame information having a frame sizefield indicating whether a storage frame is one of a small block of datain memory or a large block of data in memory. If the indicated storageframe is a small block of data, clearing the small block of data. Theclearing operation setting the bytes of the small block of data to zero.If the indicated storage frame is a large block of data, obtaining froma second general register an operand address of an initial first blockof data of the large block of data. The large block of data having afirst plurality of first blocks of data. If the indicated storage frameis a large block of data, clearing a second plurality of the firstplurality of first blocks of data, wherein the clearing sets the bytesof the second plurality of first blocks of data to zero.

In another embodiment, a block size indicator is obtained from a fieldof the machine instruction. Based on the block size indicator,determining the addressed operand is either the large block of data orthe small block of data, wherein the small block of data is the samesize as the first block of data. For the large block of data, saving anaddress of a next block of data in the second general register, the nextblock of data being a block of data following the first plurality offirst blocks of data. For the large block of data, the address of thenext block of data being determined by either encountering a boundary ofthe large block of data or responding to a program interruption event.

Perform Frame Management Function (PFMF)—Storage Keys

One embodiment of the logic of the Perform Frame Management Function inwhich the associated storage keys are set according to the instructionis described with reference to FIG. 9B.

Generally, this embodiment involves, at 950, first obtaining a machineinstruction defined for the machine architecture, which includes anopcode for a frame management instruction. The machine instruction has afirst field identifying a first general register and a second fieldidentifying a second general register. At, 952, obtaining from thesecond general register the address of an initial first block of datawithin a large block of data in main storage or memory. At 954, framemanagement information is obtained from the first general register. Theframe management information includes a key field. The key field havingfirst access protection bits. At 956, for a large block of data, asecond operand address of an initial first block of data of the largeblock of data is obtained from the second general register. The largeblock of data having a first plurality of first blocks of data. Each ofthe first plurality of first blocks having a corresponding storage keyof a first plurality of storage keys. Each storage key having storageaccess protection bits. Then, for the large block of data, the accessprotection bits of the key field are set into the storage accessprotection bits of each storage key of a second plurality of the storagekeys.

For the small block of data, an operand address of the small block ofdata is obtained from the second general register. The small block ofdata having a corresponding storage key. The storage key having storageaccess protection bits. For the small block of data, the accessprotection bits of the key field are set into the storage accessprotection bits of the storage key.

In another embodiment, a frame management field is obtained from thefirst general register. The field has a key field with a plurality ofaccess-protection bits and a block size indication. If the block sizeindication indicates a large block, a second operand address of a firstblock of data within a large block of data is obtained from a secondgeneral register The large block of data having a plurality of blocks ofdata, wherein each of the plurality of blocks of data is associated witha corresponding storage key with a plurality of storage keyaccess-protection bits. If the block size indication indicates a largeblock, the storage key access-protection bits of each correspondingstorage key of each of the plurality of small blocks of data are setwith the access-protection bits of the key field.

Alternatively, if the block size indication indicates a small block,obtaining from the second general register an operand address of thesmall block of data. The small block of data having a correspondingstorage key having storage access-protection bits. If the block sizeindication indicates a small block, the access-protection bits of thekey field are set into the storage access-protection bits of the storagekey.

In another embodiment, the block size indication is obtained from anyfield of the machine instruction, or a field of the first generalregister. Based on the block size indication, determining whether theaddressed operand is one of the large block of data or the small blockof data. The small block of data being a same size as the first block ofdata. The operand address being one of an absolute address of the largeblock of data or the real address of the small block of data. The realaddress further being subject to prefixing to determine the absoluteaddress. Further, for the large block of data, saving an address of anext block of data in the second general register, the next block ofdata being a block of data following the first plurality of first blocksof data. Additionally, for the large block of data, determining theaddress of the next block of data by any one of encountering a boundaryof the large block of data or responding to a program interruptionevent.

In yet another embodiment, the frame management field further having areference control field and a change control field. The key fieldfurther having a fetch-protection bit, a change bit and a reference bit.The storage key further having a storage fetch-protection bit, a storagereference bit and a storage change bit. For the large block of data, ifthe reference control field and the change control field are notenabled, setting the fetch-protection bit, the reference bit and thechange bit of the key field into corresponding storage fetch-protectionbits, storage reference bits and storage change bits of each storage keyof the second plurality of the storage keys. If either the referencecontrol field or the change control field are enabled and either theaccess-protection bits of the key field are not equal to the storageaccess-protection bits of the storage key or the protection bit is notequal to the storage protection bit, setting the fetch-protection bit,the reference bit and the change bit of the key field into thecorresponding storage fetch-protection bits, storage reference bits andstorage change bits of each storage key of the second plurality of thestorage keys. If the reference control field and the change controlfield are enabled, the access-protection bits of the key field beingequal to the storage access-protection bits of the storage key, and thefetch-protection bit being equal to the storage fetch-protection bit,not changing the storage fetch-protection bits, storage reference bitsand storage change bits of each storage key of the second plurality ofthe storage keys.

PFMF Optimizations for Pageable Guests

In accordance with an aspect of the present invention, the guest memorythat is the target of a PFMF instruction may or may not be backed byhost memory. If the guest memory is backed by host memory, then theinstruction performs as above. However, if the guest memory is notbacked by host memory, then certain optimizations can be performedabsent host intervention. For example, if the PFMF instruction is toperform a clear operation (e.g., logically set a guest frame to zero orzero the frame), the operation is performed without bringing in theprior host page contents from auxiliary storage to host memory. There isno reason to bring in the contents just to overlay the contents withzeros. Furthermore, the copy of the host page on auxiliary storage is nolonger needed and may be discarded. Further, if the PFMF instruction isto perform a set key operation, then that operation is optimized to usethe key stored in a host control block (i.e., the PGSTE). That is, thereis a field in the PGSTE to hold the storage key value when the frame isnot backed in host memory.

An overview of the optimizations realized when a clear operation isbeing performed is described with reference to FIG. 10. In one example,the clear operation is specified in a PFMF instruction. However, inanother embodiment, this is not so. The clear operation is specified byitself or in another instruction, etc. In the example herein, the guestframe is a host page and the page is 4K-bytes. However, in otherexamples, the guest frame may be other than a host page and/or the hostpage may be other than 4K-bytes. For example, the guest frame may be alarge frame backed by multiple small host pages, or the guest frame maybe a small frame backed by a portion of a large (for example, 1M-byte)host page. FIG. 10 illustrates a clear operation as it applies to asingle host page. If multiple host pages are involved, then the processof FIG. 10 is effectively repeated for each host page. FIG. 11, to bedescribed later, illustrates this case in detail.

Referring to FIG. 10, initially, a determination is made as to whetherthe host page containing the guest frame that is the target of the clearoperation is resident in host memory, INQUIRY 1000. If it is resident(e.g., if invalid bit 604 is off in host PTE 600), then the page iscleared, as if it was a native instruction performing the clearoperation, STEP 1002. However, if the page is not resident (i.e., theguest frame is not backed by host memory (e.g., the invalid bit is on)),then a further determination is made as to the intended usage of thepage, INQUIRY 1004. If a usage intention (UI) indicator is set to, forinstance, zero indicating that the program is not intending to use thisguest frame in the immediate future, then a host frame is not committedat this time. Instead, the following actions are taken absent hostintervention: the host page is marked logically zero by, for example,setting the Z indicator in the page status table entry corresponding tothis page, STEP 1006; the host page containing this guest frame is leftnon-resident, STEP 1008; and the guest absolute address (as an example)of the page is placed onto a backing reclaim log, so that the host knowsthat it can reclaim the backing auxiliary storage which holds the priorcontents, STEP 1010.

Returning to INQUIRY 1004, if the usage indicator is, for instance, setto one indicating that the guest frame is about to be referenced (again,as defined by the program), then the host page mapping the guest frameis made resident without host intervention through, for instance, anHPMA resolve operation. For example, a host frame is pulled off of alist of cleared available frames, STEP 1012, and is attached to the hostpage, STEP 1014. The list of cleared available frames includes one ormore pre-cleared host frames. By attaching the cleared host frame to thehost page, the guest frame mapped in that host page is backed (made hostresident) by host memory already cleared. Further details on HPMAresolve are described below with reference to FIG. 13.

Further details regarding processing a PFMF instruction (or one or moreoperations of the PFMF instruction) in a pageable guest environment,such that the performance of the operations is optimized, are describedwith reference to FIGS. 11A-11E. In particular, FIGS. 11A-11E depict oneembodiment of the logic to process a Perform Frame Management Functionby a pageable guest, in accordance with one or more aspects of thepresent invention. Although this logic describes aspects of executingthe PFMF instruction by a pageable guest, one or more aspects of thepresent invention are equally applicable to the clear frame and/or setkey operations of the instruction separate and apart from the PFMFinstruction. Further, for one or more aspects of the present invention,it is not necessary to perform both the clear frame and set keyfunctions. Aspects of the invention can apply to just one or the otheror both.

Referring to FIG. 11 a, initially various housekeeping details arehandled, in this embodiment. In other embodiments, however, one or moreof these details can be ignored or eliminated. A determination is madeas to whether the guest is in the z/Architecture® mode and whether theenhanced Dynamic Address Translation (DAT) facility (EDAT) is installed,INQUIRY 1100. If either the guest is not in z/Architecture® mode or ifEDAT is not installed, then a guest operation exception is presented,STEP 1102. In this example, PFMF is available under the enhanced DATfacility of z/Architecture®. However, this may be different for otherembodiments. One example of EDAT is described in U.S. Publication No.2009/0187724A1, entitled, “Dynamic Address Translation with FrameManagement,” Greiner et al., (IBM Docket No.: POU920070313US1),published Jul. 23, 2009, which is hereby incorporated herein byreference in its entirety.

Should EDAT be installed under the z/Architecture®, then a furtherdetermination is made as to whether the host has enabled interpretiveexecution of PFMF for this guest, INQUIRY 1104. As one example, this isindicated by a bit in a control block, known as a state description,that is created by the host to represent the state of a virtual machineand is used by the firmware in interpretive execution.

If the host has not enabled interpretive execution, then an instructioninterception is presented to the host, STEP 1106. This allows the hostto trap the instruction, if desired, for debugging and special-casesimulation.

However, if interpretive execution is enabled, then a check is made tosee if the guest program status word (PSW) specifies problem state,INQUIRY 1108. If so, then a guest privileged operation exception ispresented, STEP 1110. In this embodiment, PFMF is a privilegedinstruction for use by operating systems, not application programs. Thismay be different, however, in other embodiments.

Should the PSW not specify problem state, then a determination is madeas to whether the reserved bits are non-zero in general register R1(i.e., the general register whose number appears as the R1 operand inthe instruction), or if the frame size code (FSC) in R1 has a valueother than binary 000 or 001, INQUIRY 1112. If the reserved bits arenon-zero or the frame size code has a value other than binary 000 or001, then a guest specification exception is presented, STEP 1114.

Otherwise, the guest absolute address is obtained, STEP 1116. In oneexample, to obtain the guest absolute address, the frame size code inregister R1 is examined. If FSC is binary 000, then the operand is a4K-byte guest frame specified by the guest real address in R2. Guestprefixing is applied to derive the target guest absolute address. If FSCis binary 001, then the operand is a 1M-byte frame specified by theguest absolute address in R2. This address is used as is.

Thereafter, an interval completion indicator in the state description isset and a checkpoint synchronization operation is performed, STEP 1118.This is done so that if a machine failure occurs during subsequentprocessing, the host can recognize that host data structures may havebeen corrupted.

For each 4K-byte block of guest storage designated by the operand, i.e.,for a single 4K-byte frame if FSC is binary 000 or for each of the4K-byte blocks beginning with the operand address and continuing to thenext 1M-byte boundary if FSC is binary 001, the following steps areperformed, STEP 1120:

-   -   The guest absolute address of the frame is treated as a host        virtual address, and host dynamic address translation is        performed on that address to locate the leaf host DAT-table        entry (either a valid format-1 segment table entry designating a        large host frame or a page table entry indicating a small frame        size), STEP 1122. Examples of DAT are described in        “z/Architecture Principles of Operation,” IBM® Publication No.        SA22-7832-05, April, 2007; and U.S. Publication No.        2009/0187724A1, entitled, “Dynamic Address Translation with        Frame Management,” Greiner et al., (IBM Docket No.:        POU920070313US1), published Jul. 23, 2009, each of which is        hereby incorporated herein by reference in its entirety.    -   Various error processing associated with DAT is described with        reference to FIG. 12. For instance, if a host ASCE-type, region        translation or segment translation exception condition is        encountered preventing the leaf host DAT table entry for the        translation from being located, INQUIRY 1200, then checkpoint        synchronization is performed and an interval completion        indicator is turned off, STEP 1201. This indicates to the host        that updates to the host data structures are complete, and the        structures are in a consistent state should a machine failure        occur. Additionally, the remainder of instruction execution is        suppressed and an instruction interception is presented to the        host, STEP 1202. This allows the host to resolve the fault        condition and at the same time, optimize the handling of PFMF        for the case where there is no prior backing content and no host        page table or page status table is available.    -   Otherwise, in one embodiment, if FSC is binary 001, specifying a        large (e.g., 1 MB) guest frame, and storage is mapped by small        host pages (format control (FC) in host segment table entry is        zero), INQUIRY 1204, then checkpoint synchronization is        performed and an interval completion indicator is turned off,        STEP 1201. Further, an instruction interception is presented to        the host, STEP 1202. This allows the host to optimize handling        by discarding any prior small-frame backing and backing with a        large host frame, if desired. In other embodiments, this test        may be omitted or made conditional on, for example, a control        bit in the state description.    -   If the case tested at INQUIRY 1204 does not apply (or if, in        another embodiment, the test is bypassed), a determination is        next made as to whether host DAT protection is indicated in any        of the host RTEs, STE or PTE used in the translation, and either        the clear frame (CF) or the set key (SK) operation is requested        in the R1 operand, INQUIRY 1206. For the PFMF instruction, DAT        protection is recognized even if the leaf host table entry is        invalid, to prevent handling which would alter the content of        the host page or storage key in the non-host-resident case. If        DAT protection is indicated, then checkpoint synchronization is        performed and an interval completion indicator is turned off,        STEP 1207. Additionally, a protection exception is presented,        STEP 1208.    -   Otherwise, if any other access exception is encountered, such as        an invalid host address, INQUIRY 1210, then checkpoint        synchronization is performed and an interval completion        indicator is turned off, STEP 1211. Additionally, this access        exception is presented, as usual, STEP 1212.    -   Returning to FIG. 11A, subsequent to performing host dynamic        address translation to locate the leaf host DAT table entry, a        determination is made as to whether the host ASCE was a        real-space designation (RSD), for which there are no table        entries, or whether the leaf entry is a format-1 STE, INQUIRY        1124. In either of these cases, the guest frame is resident in        host memory, so optimizations are not needed. A clear frame        and/or set key function, as requested in the CF and SK        indicators of the R1 operand, is/are performed directly into the        designated 4K-byte block within the host memory containing the        guest frame, STEP 1126. In the case of a host RSD, this is the        4K-byte block at the host real address matching the guest        absolute address. In the case of a format-1 STE, this is the        large frame designated by the segment frame absolute address        field of this STE. Processing then continues by advancing to the        next 4K-byte block, if any, STEP 1166 (FIG. 11D), described        below.    -   On the other hand, if the host ASCE is not an RSD and the leaf        entry is not a format-1 STE, INQUIRY 1124, then the leaf entry        is a PTE and processing continues, as described below.    -   The PGSTE associated with the host PTE is located, STEP 1128,        for example, by adding a fixed offset to the PTE address to        obtain the PGSTE address. Next, a determination is made as to        whether the PGSTE is at an invalid address, INQUIRY 1130. If so,        then a checkpoint-synchronization operation is performed and the        interval completion indicator is turned off, STEP 1131. Further,        instruction execution is terminated and a validity interception        is presented to the host, STEP 1132. This signifies an error in        the host data structures.    -   If the PGSTE address is valid, INQUIRY 1130, an attempt is made        to obtain the page control lock (PCL) in the host PGSTE        corresponding to the host PTE, STEP 1133. In one example, this        attempt is made by changing the PCL bit from zero to one using        an interlocked update. If this operation fails because the PCL        bit is already one, INQUIRY 1134, then a checkpoint        synchronization operation is performed and the interval        completion indicator is turned off, STEP 1135. Moreover, the        remainder of the instruction is suppressed and an instruction        interception is presented to the host, STEP 1136.    -   Should the PCL be obtained successfully, INQUIRY 1134, then the        host PTE is fetched and examined, STEP 1138. If the invalid bit        in the PTE is off, signifying that the page is resident in host        storage, INQUIRY 1140, then a clear frame and/or set key        function, as requested in the CF and SK indicators of the R1        operand, is/are performed directly into the designated host        small (4K-byte) frame designated by the page frame real address        field of this PTE, STEP 1142. Processing then proceeds to        INQUIRY 1160 (FIG. 11D), as described below.    -   Returning to INQUIRY 1140, if the invalid bit is on, and a clear        frame function is requested (i.e., the CF indicator is on),        INQUIRY 1144 (FIG. 11C), then processing of the clear frame        operation proceeds according to one of the following scenarios:        -   (1) If the usage indicator in the R1 operand is on (e.g.,            set to 1) specifying that the guest has indicated that it            intends to make use of the frame contents in the near            future, INQUIRY 1146, then a Host Page Management Assist            (HPMA) resolve operation is performed to make the host page            resident with zero contents, STEP 1148.            -   One example of the HPMA operation is described in U.S.                Publication No. 2005/0268071A1, entitled “Facilitating                Management of Storage of a Pageable Mode Virtual                Environment Absent Intervention of a Host of the                Environment,” Blandy et al., published Dec. 1, 2005,                which is hereby incorporated herein by reference in its                entirety. Moreover, one example of HPMA resolve                processing, in accordance with an aspect of the present                invention, is described further below with reference to                FIGS. 13A-13B.            -   As described with reference to FIGS. 13A-13B, if HPMA                resolve processing fails, then processing of the guest                PFMF instruction stops at this point, with an                instruction interception or validity interception                presented to the host. Otherwise, processing continues                at INQUIRY 1160 (FIG. 11D), described below.        -   (2) If the usage indicator in the R1 operand is off (e.g.,            set to 0), INQUIRY 1146, and the logically zero (Z) bit in            the PGSTE is off, INQUIRY 1150, then the guest absolute            (i.e., host virtual) address of this 4K-byte block is            appended to the CMM (Collaborative Memory Management)            backing reclaim log (CBRL) designated by the state            description, STEP 1152. One example of the CBRL and            associated processing is described in U.S. Patent            Application Publication No. US 2007/0016904 A1, entitled,            “Facilitating Processing Within Computing Environment            Supporting Pageable Guests,” Adlung et al., published Jan.            18, 2007, which is hereby incorporated herein by reference            in its entirety. Moreover, one example of appending to the            CBRL, in accordance with an aspect of the present invention,            is described further below with reference to FIG. 14.            -   As described with reference to FIG. 14, if appending to                the CBRL fails, then processing of the guest PFMF                instruction stops at this point, with an instruction                interception or validity interception presented to the                host. Otherwise, the Z bit, which causes subsequent host                or firmware actions to treat the host page contents as                zeros and allows HPMA resolve to provide a host frame of                zeros when the guest later references this storage, is                set, STEP 1153, and processing continues at INQUIRY                1154, as described below.        -   (3) If the usage indicator is off, INQUIRY 1146, and the Z            bit is already on, INQUIRY 1150, then no action is needed to            clear the contents. The page is already marked as logically            zero. Processing continues at INQUIRY 1154, as described            below.    -   If the invalid bit is on, INQUIRY 1140 (FIG. 11B) and the set        key function is requested (i.e., the SK bit is on), INQUIRY 1154        (FIG. 11C), then the access control and fetch protection (ACC,F)        values 826, 828 from the R1 operand are placed into the ACC and        FP fields 702, 704, respectively, of the PGSTE, STEP 1156, and        guest reference and guest change fields 712, 714 in the PGSTE        are set based on the R and C values 830, 832 specified in the R1        operand, STEP 1158. This records the requested key value with        the host page, so that it can be interrogated in the PGSTE,        while the host page is not host resident, and set into the host        frame when the page is later made resident.    -   If either the clear frame or the set key function was requested,        INQUIRY 1160 (FIG. 11D), then the CMM block usage state in the        PGSTE is set to stable (e.g., binary 00), STEP 1162. This is the        appropriate setting for a guest frame being put into use.        Handling this here takes advantage of the PCL already being held        and avoids a separate instruction (e.g., ESSA instruction) by        the guest.    -   Thereafter, or if neither clear nor set was requested, the PCL        is released, STEP 1164, by, for instance, setting the PCL bit in        the PGSTE to zero, and processing for this 4K-byte block is        complete.    -   Next, a determination is made as to whether a large (e.g.,        1M-byte) guest frame size was specified (e.g., FSC=001), INQUIRY        1166. If so, then 4K is added to the operand address in R2, STEP        1168, and a determination is made as to whether this addition        resulted in an address beyond the end of the guest frame        specified as the operand, INQUIRY 1170. That is, if the addition        results in a carry into the next megabyte address, then the loop        to process each 4K-byte block in the 1M-byte guest frame, which        began at STEP 1120, is exited. If the end of the guest frame has        not yet been reached, then a determination is next made as to        whether an asynchronous interruption is pending, INQUIRY 1172.        If an asynchronous interruption is pending, the unit of        operation is nullified, STEP 1174. For instance, the instruction        address in the guest PSW is backed up to point to this PFMF        instruction. Thereafter, the loop is exited. This allows the        interruption to be taken, after which the PFMF instruction can        be re-executed and will resume with the next 4K-byte block to be        processed.    -   If there is no asynchronous interruption pending, INQUIRY 1172,        then processing returns to STEP 1120 and the logic is repeated        with the next 4K-byte block.

After exiting the loop, a checkpoint synchronization operation isperformed, STEP 1180 (FIG. 11E), and the interval completion indicatorin the state description is turned off, STEP 1182. This indicates to thehost that updates to the host data structures are complete, and they arein a consistent state should a machine failure occur. Processing of theguest PFMF instruction is now complete.

In the above description, reference is made to HPMA resolve processing.Further details regarding one embodiment of this processing aredescribed with reference to FIGS. 13A-13B. In this embodiment, HPMAresolve is used to assign host frames to host pages. Initially, theanchor of the available frame descriptor list (e.g., at location 210 hexin the host prefix area) is examined, STEP 1300. If the anchor is zeroindicating that the list is empty, INQUIRY 1302, then the PCL isreleased, a checkpoint synchronization operation is performed, and theinterval completion indicator is turned off, STEP 1304. In one example,the PCL is released by setting the PCL bit in the PGSTE to zero.Additionally, the remainder of the instruction processing is suppressed,STEP 1306, and an instruction interception is presented to the host,STEP 1308. This allows the host to complete the execution of the guestPFMF, as well as to replenish the available frame list for future use.

Returning to INQUIRY 1302, if the anchor is not zero, the host frameaddress is extracted from the frame descriptor at the front of the list(i.e., the one to which the anchor points), STEP 1310. This is anaddress of a frame available for use whose contents the host has alreadycleared to zeros.

If the set key function is requested (i.e., the SK indicator is on),INQUIRY 1312, then the access control and fetch protection field in thestorage key of this host frame is set based on the key value specifiedin the R1 operand, STEP 1314. Further, the reference and changeindicators in the storage key are set to zeros, STEP 1316, and the guestreference and guest change indicators in the PGSTE are set to thecorresponding values specified in the R1 operand, STEP 1318.

Returning to INQUIRY 1312, if the set key function is not requested, thestorage key of the frame is set to the value already present in theaccess control and fetch protection (ACC,FP) fields of the PGSTE, STEP1320, and the reference and change bits in the storage key are set tozeros, STEP 1322. This ensures that the guest view of the storage keywill have the same value when the block is host resident that it hadbefore the block was made resident.

In addition to setting the change bits (STEP 1318 or 1322), the firstframe descriptor is dequeued from the available frame descriptor list bycopying its forward pointer to the anchor, STEP 1324 (FIG. 13B).

Moreover, the assignment of the host frame to the host page is recordedby, for instance, copying the current PTE contents into the PTE copyfield of the frame descriptor and storing the PTE address into the PTEaddress field of the frame descriptor, STEP 1326. This allows the hostto update its data structures at a later point and to locate informationdescribing backing storage for the host page that can be reclaimed.

The page content replacement indicator is set in the PGSTE to indicatethat this is a resolve operation with content replacement, STEP 1328.This allows the host to keep statistics on such operations. Further, thepage on processed list indicator is set in the PGSTE to indicate thatthe frame assigned to this page may be found on the processed framedescriptor list, STEP 1330.

The PTE is updated to remove the page invalid condition, STEP 1332. Inone example, this is accomplished by storing revised contents includinga page frame real address equal to the extracted frame address, aninvalid bit with a value of zero, and remaining fields, including theDAT protection bit, copied unchanged from the old PTE contents.

The frame descriptor is added to the front of the processed framedescriptor list, anchored at, for example, location 210 hex in the hostprefix area, STEP 1334. This is accomplished by, for instance, storingthe prior contents of this anchor into the forward pointer in the framedescriptor, and then performing an interlocked update to replace theprior anchor contents with the address of the frame descriptor. If theinterlocked update fails because the anchor contents have changed duringthis step, then the step is repeated until it succeeds. This concludesthe processing associated with the HPMA resolve operation.

Reference is also made in the above description to processing associatedwith appending to the CMM backing reclaim log (CBRL). One embodiment ofthe logic associated with this processing is described with reference toFIG. 14. Initially, a determination is made as to whether the CBRLorigin in the state description is zero, INQUIRY 1400. If the CBRLorigin is zero signifying that there is no CBRL, then the PCL isreleased, a checkpoint synchronization operation is performed, and theinterval completion indicator is turned off, STEP 1402. Further, theremainder of the PFMF instruction execution is suppressed, STEP 1404,and an instruction interception is presented to the host, STEP 1406.This allows the host to allocate memory for a CBRL or to handle the PFMFin simulation.

If the CBRL origin is not zero, but it designates an invalid address,INQUIRY 1408, the PCL is released, a checkpoint synchronizationoperation is performed, and the interval completion indicator is turnedoff, STEP 1410. Moreover, the remainder of the instruction execution issuppressed, STEP 1412, and a validity interception is presented to thehost, STEP 1414. This signifies an error in the host data structures.

Should the CBRL origin designate a valid address, a determination ismade as to whether the CBRL next entry offset (NEO) is a given value(e.g., hex FF8), signifying that the CBRL is full, INQUIRY 1416. If so,the PCL is released, a checkpoint synchronization operation isperformed, and the interval completion indicator is turned off, STEP1402. Additionally, the remainder of instruction execution issuppressed, STEP 1404, and an instruction interception is presented tothe host, STEP 1406. This allows the host to process the CBRL and handlethe PFMF in simulation.

Otherwise, if the CBRL is not full, the guest absolute block address isstored into the CBRL at the location specified by the NEO, STEP 1418.Further, the NEO is incremented by the length of the address stored, forinstance, eight bytes, STEP 1420. This adds the address to the log. Thehost can subsequently process the log in order to release backingstorage resources including the prior page contents, which are no longerneeded because the PFMF clear frame operation rendered the contentslogically zero. This concludes one example of the CBRL processing.

Described in detail above are optimizations provided for a clear and/orset operation when issued by a pageable guest. As one example,optimizations are provided for operations (e.g., clear frame and/or setkey) of a Perform Frame Management Function.

One or more aspects of the present invention can be included in anarticle of manufacture (e.g., one or more computer program products)having, for instance, computer usable media. The media has therein, forinstance, computer readable program code means or logic (e.g.,instructions, code, commands, etc.) to provide and facilitate thecapabilities of the present invention. The article of manufacture can beincluded as a part of a computer system or sold separately.

One example of an article of manufacture or a computer program productincorporating one or more aspects of the present invention is describedwith reference to FIG. 15. A computer program product 1500 includes, forinstance, one or more computer usable media 1502 to store computerreadable program code means or logic 1504 thereon to provide andfacilitate one or more aspects of the present invention. The medium canbe an electronic, magnetic, optical, electromagnetic, infrared, orsemiconductor system (or apparatus or device) or a propagation medium.Examples of a computer readable medium include a semiconductor or solidstate memory, magnetic tape, a removable computer diskette, a randomaccess memory (RAM), a read-only memory (ROM), a rigid magnetic disk andan optical disk. Examples of optical disks include compact disk-readonly memory (CD-ROM), compact disk-read/write (CD-R/W) and DVD.

A sequence of program instructions or a logical assembly of one or moreinterrelated modules defined by one or more computer readable programcode means or logic direct the performance of one or more aspects of thepresent invention.

Advantageously, various optimizations are provided, in accordance withone or more aspects of the present invention. For example, a guestrequest to clear a 4K-byte frame which is not host-resident may besatisfied without a transition to the host, either by binding the hostpage mapping the guest frame to a cleared host frame from a listpreviously supplied by the host (HPMA), or by marking the host pagelogically zero and appending its address to a CMM-backing-reclaim log(CMM). Which of these actions is taken depends on a usage-intentindicator specified by the guest. Similarly, setting of the guest key isoptimized for a guest frame that is not host resident by placing the keyvalue directly into the host PGSTE.

Moreover, the following advantages are provided, in accordance with oneor more aspects of the present invention:

-   -   Host page faults and page reads of old guest memory contents        which are about to be cleared are avoided.    -   If the guest memory to be cleared is not host resident, but is        expected to be referenced soon (per the UI bit), the guest        memory addresses are bound to pre-cleared host memory frames        from an available list supplied in advance by the host, thus        avoiding the overhead of clearing in-line and avoiding a context        switch to and from the host program.    -   If the guest memory to be cleared is not host resident and is        not expected to be referenced soon, the host page containing the        guest memory is left invalid (i.e., non-resident) and is marked        logically zero. This avoids context switches and clearing at the        time of the PFMF. In this case, the host page is also added to a        backing-reclaim log, that is, a list of host pages whose        auxiliary backing storage can be reclaimed. This allows the host        to reuse the backing resources more timely.    -   Following the logically zero operation above, the host page is        in a state wherein a subsequent reference to the guest frame can        also be resolved without context switches, using the pre-cleared        available list described above.    -   If PFMF requests setting of the storage key for a guest memory        mapped by a host page which is and remains not host-resident        (i.e., if the page was not resident, and either no clearing was        requested, or UI was zero), the key value is placed into the        host page-status-table entry for the page, as is done when Set        Storage Key Extended (SSKE) is performed on a host-non-resident        page. This allows the proper key value to be retained without        requiring a physical storage key, which is only available to a        resident page.    -   If set key is performed in the same operation as a clear frame,        host serialization can be obtained once instead of twice.    -   A PFMF operation which requests either frame clearing or key        setting also changes the collaborative-memory-management state        of the guest frame to stable. This is the desired effect when a        guest memory is being redeployed, and eliminates the need for        the guest to issue separate ESSA instruction(s).    -   In one embodiment, the above enhancements operate on host page        metadata which is available for 4 KB host pages. When the guest        invokes PFMF specifying a 1 MB frame, the operation is instead        intercepted to the host, which can then re-back the guest frame        with a host 1 MB frame if desired. In this case, the host need        not page in and copy the old frame contents, but can simply        clear a new 1 MB host frame (or use one from a pre-cleared        list). Alternatively, in another embodiment, processing of a        guest 1 MB frame mapped by host 4 KB pages is performed on the        individual pages, avoiding the interception for example, for a        host that does not employ 1 MB host frames.    -   If a higher-level invalid host table entry (a region- or        segment-table entry) prevents access to the host page table and        page status table, PFMF is again intercepted to the host, rather        than presenting an ordinary region- or segment-translation        exception. This allows the host to use the page size specified        on PFMF as a hint, and to allocate host backing memory of the        same size if desired, or to delay allocation of backing memory        if the usage intent is not indicated.

In a further embodiment, metadata at the host segment level (e.g., for 1MB units of host memory) is provided. In this case, PFMF interpretiveexecution could take advantage of these to avoid interception when guestframe size and host page size coincide.

Commercial Implementation

Although the z/Architecture® by IBM® is mentioned herein, one or moreaspects of the present invention are equally applicable to other machinearchitectures and/or computing environments employing pageable entitiesor similar constructs.

Commercial implementations of the PFMF instruction, facilities, andother formats, instructions, and attributes disclosed herein can beimplemented either in hardware or by programmers, such as operatingsystem programmers, writing in, for example, assembly language. Suchprogramming instructions may be stored on a storage medium intended tobe executed natively in a computing environment, such as thez/Architecture® IBM® Server, or alternatively in machines executingother architectures. The instructions can be emulated in existing and infuture IBM® servers and on other machines or mainframes. They can beexecuted in machines where generally execution is in an emulation mode.

In emulation mode, the specific instruction being emulated is decoded,and a subroutine is built to implement the individual instruction, as ina subroutine or driver, or some other technique is used for providing adriver for the specific hardware, as is within the skill of those in theart after understanding the description hereof. Various software andhardware emulation techniques are described in numerous United Statespatents including: U.S. Pat. Nos. 5,551,013, 5,574,873, 5,790,825,6,009,261, 6,308,255, and 6,463,582, each of which is herebyincorporated herein by reference in its entirety. Many other teachingsfurther illustrate a variety of ways to achieve emulation of aninstruction format architected for a target machine.

Other Variations And Architectures

While various examples and embodiments are described herein, these areonly examples, and many variations are included within the scope of thepresent invention. For example, the computing environment describedherein is only one example. Many other environments, including othertypes of communications environments, may include one or more aspects ofthe present invention. For instance, different types of processors,guests and/or hosts may be employed. Moreover, other types ofarchitectures can employ one or more aspects of the present invention.

Aspects of the invention are beneficial to many types of environments,including environments that have a plurality of zones, andnon-partitioned environments. Further, there may be no central processorcomplexes, but yet, multiple processors coupled together. Variousaspects hereof are applicable to single processor environments.

Further, in the examples of the data structures and flows providedherein, the creation and/or use of different fields may include manyvariations, such as a different number of bits; bits in a differentorder; more, less or different bits than described herein; more, less ordifferent fields; fields in a differing order; different sizes offields; etc. Again, these fields were only provided as an example, andmany variations may be included. Further, indicators and/or controlsdescribed herein may be of many different forms. For instance, they maybe represented in a manner other than by bits. Additionally, althoughthe term address is used herein, any designation may be used.

As used herein, the term “page” is used to refer to a fixed-size orpredefined-size area of virtual storage (i.e., virtual memory). As oneexample, a host page is an area of host virtual storage. The size of thepage can vary, although in the examples provided herein, a page is 4Kbytes. Further, a “frame” is used to refer to a fixed-size or predefinedsize area of real or absolute storage (i.e., memory). As examples, ahost frame is an area of host real or absolute storage, and a guestframe is an area of guest real or absolute storage. In the case of apageable guest, this guest real or absolute storage is mapped by hostvirtual storage. As is common, pages of host virtual storage are backedby frames of host real or absolute storage, as needed. The size of theframe can vary, although in the examples provided herein, a frame is4K-bytes or 1M-bytes. However, in other embodiments, there may bedifferent sizes of pages, frames, segments, regions, blocks of storage,etc. Moreover, in other architectures, the terms “page” and “segment”may be used interchangeably or the term “page” may be used to apply tomultiple size units of virtual storage. The term “obtaining”, such asobtaining an instruction, includes, but is not limited to, fetching,having, receiving, being provided, creating, forming, issuing, etc. Aninstruction can reference other registers or can reference other thanregisters, such as operands, fields, locations, etc. Many otheralternatives to the above are possible. Further, although terms, such aslists, tables, etc. are used herein, any types of data structures may beused. For instance, a table can include other data structures as well.Again, those mentioned herein are just examples.

Further, a data processing system suitable for storing and/or executingprogram code is usable that includes at least one processor coupleddirectly or indirectly to memory elements through a system bus. Thememory elements include, for instance, local memory employed duringactual execution of the program code, bulk storage, and cache memorywhich provide temporary storage of at least some program code in orderto reduce the number of times code must be retrieved from bulk storageduring execution.

Input/Output or I/O devices (including, but not limited to, keyboards,displays, pointing devices, DASD, tape, CDs, DVDs, thumb drives andother memory media, etc.) can be coupled to the system either directlyor through intervening I/O controllers. Network adapters may also becoupled to the system to enable the data processing system to becomecoupled to other data processing systems or remote printers or storagedevices through intervening private or public networks. Modems, cablemodems, and Ethernet cards are just a few of the available types ofnetwork adapters.

The capabilities of one or more aspects of the present invention can beimplemented in software, firmware, hardware, or some combinationthereof. At least one program storage device readable by a machineembodying at least one program of instructions executable by the machineto perform the capabilities of the present invention can be provided.

The flow diagrams depicted herein are just examples. There may be manyvariations to these diagrams or the steps (or operations) describedtherein without departing from the spirit of the invention. Forinstance, the steps may be performed in a differing order, or steps maybe added, deleted, or modified. All of these variations are considered apart of the claimed invention.

Although embodiments have been depicted and described in detail herein,it will be apparent to those skilled in the relevant art that variousmodifications, additions, substitutions and the like can be made withoutdeparting from the spirit of the invention and these are thereforeconsidered to be within the scope of the invention as defined in thefollowing claims.

What is claimed is:
 1. A computer system to facilitate management offrames within a computing environment, said computer system comprising:a memory; and a processor in communication with the memory, wherein thecomputer system is configured to perform a method, the methodcomprising: based on initiating execution of an instruction issued by apageable guest executing within the computing environment and specifyinga clear operation, determining a usage intent characteristic for a guestframe to be cleared, the guest frame not resident in host memory; andperforming the clear operation absent host intervention, the performingbeing based on the determined usage intent characteristic.
 2. Thecomputer system of claim 1, wherein the determining the usage intentcharacteristic comprises employing an indicator of the instruction. 3.The computer system of claim 1, wherein the instruction comprises aperform frame management instruction.
 4. The computer system of claim 1,wherein the usage intent characteristic specifies whether a program hasindicated that it is likely to use the guest frame within a near future.5. The computer system of claim 4, wherein the usage intentcharacteristic specifies that the program has indicated that it islikely to use the guest frame within the near future, and wherein theperforming the clear operation comprises: obtaining a host frame from alist of cleared available frames; and attaching the obtained host frameto the guest frame to be cleared.
 6. The computer system of claim 4,wherein the usage intent characteristic specifies that the program hasindicated that it is not likely to use the guest frame within the nearfuture, and wherein the performing the clear operation comprises:marking one or more host pages that back the guest frame as logicallyzero.
 7. The computer system of claim 6, wherein the guest frame to becleared remains not host resident.
 8. The computer system of claim 1,wherein the method further comprises setting a storage key of the guestframe.
 9. The computer system of claim 8, wherein the setting thestorage key comprises including a value of the storage key in a controlblock used by a host of the computing environment, the host managing thepageable guest.
 10. The computer system of claim 9, wherein the controlblock comprises a page status table entry.
 11. A computer programproduct for facilitating management of frames within a computingenvironment, said computer program product comprising: a computerreadable storage medium readable by a processing circuit and storinginstructions for execution by the processing circuit for performing amethod comprising: based on initiating execution of an instructionissued by a pageable guest executing within the computing environmentand specifying a clear operation, determining a usage intentcharacteristic for a guest frame to be cleared, the guest frame notresident in host memory; and performing the clear operation absent hostintervention, the performing being based on the determined usage intentcharacteristic.
 12. The computer program product of claim 11, whereinthe usage intent characteristic specifies whether a program hasindicated that it is likely to use the guest frame within a near future.13. The computer program product of claim 12, wherein the usage intentcharacteristic specifies that the program has indicated that it islikely to use the guest frame within the near future, and wherein theperforming the clear operation comprises: obtaining a host frame from alist of cleared available frames; and attaching the obtained host frameto the guest frame to be cleared.
 14. The computer program product ofclaim 12, wherein the usage intent characteristic specifies that theprogram has indicated that it is not likely to use the guest framewithin the near future, and wherein the performing the clear operationcomprises: marking one or more host pages that back the guest frame aslogically zero.
 15. The computer program product of claim 11, whereinthe method further comprises setting a storage key of the guest frame.16. A method of facilitating management of frames within a computingenvironment, said method comprising: based on initiating execution of aninstruction issued by a pageable guest executing within the computingenvironment and specifying a clear operation, determining a usage intentcharacteristic for a guest frame to be cleared, the guest frame notresident in host memory; and performing the clear operation absent hostintervention, the performing being based on the determined usage intentcharacteristic.
 17. The method of claim 16, wherein the usage intentcharacteristic specifies whether a program has indicated that it islikely to use the guest frame within a near future.
 18. The method ofclaim 17, wherein the usage intent characteristic specifies that theprogram has indicated that it is likely to use the guest frame withinthe near future, and wherein the performing the clear operationcomprises: obtaining a host frame from a list of cleared availableframes; and attaching the obtained host frame to the guest frame to becleared.
 19. The method of claim 17, wherein the usage intentcharacteristic specifies that the program has indicated that it is notlikely to use the guest frame within the near future, and wherein theperforming the clear operation comprises: marking one or more host pagesthat back the guest frame as logically zero.
 20. The method of claim 16,further comprising setting a storage key of the guest frame.